Persistent Data Structures and Planar Point Location


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1 Persistent Data Structures and Planar Point Location Inge Li Gørtz
2 Persistent Data Structures Ephemeral Partial persistence Full persistence Confluent persistence V1 V1 V1 V1 V2 q ue V2 V2 V5 V2 V4 V4 r i e s V4 V7 V3 V4 V6 V8 update and query all versions V6 V7 V3 V5 V3 V3 V5 update and query last version V5 update update, query and combine all versions
3 Persistent Data Structures Ephemeral data structures. A modification destroys the old version. Persistent data structures. Modifications nondestructive. Each modification creates a new version. All versions coexists. Partially persistent data structures. Can access all versions, but only update the latest version. Fully persistent data structures. Can access and modify all versions (branching). An update operation only operates on a single version at a time (cannot combine several versions). Confluent persistent data structures. Can access and modify all versions and combine versions into new versions. Functional data structures. Cannot change any nodes, only create new ones.
4 Simple methods for making data structures persistent Structurecopying method. Create a copy of the data structure each time it is changed. Slowdown of Ω(n) time and space per update to a data structure of size n. Store a logfile of all updates. In order to access version i, first carry out i updates, starting with the initial structure, and generate version i. Overhead of Ω(i) time per access, O(1) space and time per update. Hybridmethod. Store the complete sequence of updates and additionally each kth version for a suitably chosen k. Result: Any choice of k causes blowup in either storage space or access time.
5 Overview Partial persistence. Fat node method. Node copying Full persistence. Main idea. Algorithmic applications
6 Partial Persistence Fat node method
7 Fat node method Associate set c(x) for each location in memory x. c(x)={<t,v>: x modified in version t, x has value v after construction of version t} x D(x): data structure containing c(x) Query q(t,x): Find largest version number t in t such that t t. Return value associated with t in D(x). Update (create new version m): If memory locations x1,...,xk modified to the values v1,...vk: Insert <m,vi> in D(xi).
8 Fat node method Implementation of D(x): Balanced binary search tree: query O(log c(x) ) = O(log m), m number of versions. Update: O(1) Extra space: O(1) yfast trie: query: O(loglog m) update: expected O(loglog m) Extra space: O(1)
9 Fat node method Linked data structures: each pointer field store many time value pairs. new node created by ephemeral update: create new node and mark all fields with version i. Auxiliary array keep pointer to root of each version Value 10 2
10 Fat node method example root
11 Fat node method Driscoll, Sarnak, Sleator, Tarjan, Any data structure can be made partially persistent with slowdown O(log m) for queries and O(1) for updates. The space cost is O(1) for each ephemeral memory modification. Any data structure can be made partially persistent on a RAM with slowdown O(loglog m) for queries and expected slowdown O(loglog m) for updates. The space cost is O(1) for each ephemeral memory modification
12 Partial Persistence Node copying method
13 Node copying method Linked data structure with bounded indegree p, p = O(1). Each node has p predecessor pointers + p + 1 extra fields. Auxiliary array to keep pointer to root of each version (field name, version)
14 Partially persistent balanced search trees via node copying Suffices to allow one extra pointer field in each node Each extra pointer is tagged with a version number and a field name. When ephemeral update allocates a new node: allocate a new node as well. When ephemeral update changes a pointer field: if the extra pointer is empty use it, otherwise copy the node. Try to store pointer to the new copy in its parent. If the extra pointer at the parent is occupied copy the parent and continue going up this way. Maintain array of roots indexed by timestamp. Value left 4
15 Node copying method example R L
16 Partially persistent BST with node copying Analysis: Time slowdown: access: O(1) updates: O(1) amortized Extra space: O(1) amortized O(1) for new nodes also created by ephemeral data structure O(1) amortized space for nodes created when a node is full. Proof uses potential analysis (next time).
17 Partially Persistent Data Structures Driscoll, Sarnak, Sleator, Tarjan, Any boundeddegree linked data structure can be made partially persistent with (worstcase) slowdown O(1) for queries, amortized slowdown O(1) for updates, and amortized space cost O(1) per memory modification.
18 Full Persistence Fat node method
19 Version tree Version tree: partial order. V1 Tree color problem: AddLeaf(v, c): Add leaf u as child of v, with color(u)=c. V2 V5 Lookup(v, c): Find nearest ancestor of v with color c. V3 V4 V6 V8 Fully persistent array: V7 Store(A, i, x, t): Set A[i]=x at time t ~ AddLeaf(t, i), value v = x. Access(A, i, t): Lookup value A[i] at time t ~ Lookup(t, i)
20 Version tree and version list Euler tour: L(T) = (v1, v2, v3, v7, v7, v3, v4, v4, v6, v6, v2, v5, v8, v8, v5, v1 ) Partition list for each color: V1 L(1) = (v1, v2, v3, v7, v7, v3, v4, v4 ), (v6, v6 ), (v2, v5, v8, v8, v5, v1 ) V2 V5 L(2) = (v1), (v2), (v3, v7, v7, v3 ), (v4, v4 ), (v6, v6, v2, v5, v8, v8, v5, v1 ) V3 V4 V6 V8 L(3) = (v1, v2, v3, v7, v7, v3, v4, v4, v6, v6, v2 ), (v5), (v8, v8 ), (v5, v1 ) V7 Predecessor data structure for each color to find right sublist. Maintaining order in a list problem: O(1) time.
21 Fully Persistent Data Structures Driscoll, Sarnak, Sleator, Tarjan, Any data structure can be made fully persistent with slowdown O(log m) for both queries and updates. The space cost is O(1) for each ephemeral memory modification. Any boundeddegree linked data structure can be made fully persistent with (worstcase) slowdown O(1) for queries, amortized slowdown O(1) for updates, and amortized space cost O(1) per memory modification. Dietz, Any data structure can be made fully persistent on a RAM with slowdown O(loglog m) for queries and expected slowdown O(loglog m) for updates. The space cost is O(1) for each ephemeral memory modification.
22 Algorithmic Applications
23 Planar Point Location Planar point location. Euclidean plane subdivided into polygons by n line segments that intersect only at their endpoints. Query: given a query point p determine which polygon that contains p. Measure algorithm by three parameters: Preprocessing time Query time Space
24 Planar point location: Example From slides by H. Kaplan
25 Planar point location: Example From slides by H. Kaplan 25
26 Planar Point Location Within each slab the lines are totally ordered. Search tree per slab containing the lines at the leaves with each line associate the polygon above it. Another search tree on the xcoordinates of the vertical lines. query find appropriate slab search the search tree of the slab to find the polygon
27 Planar Point Location One search tree for each slab: Query time: O(log n) Space: Ω(n 2 ) Total # lines O(n), and number of lines in each slab is O(n).
28 Planar point location: Improve space bound Key observation: The lists of the lines in adjacent slabs are very similar. Create the search tree for the first slab. Obtain the next one by deleting the lines that end at the corresponding vertex and adding the lines that start at that vertex. Number of insertions/deletions? 2n Use partially persistent search tree. x axis is time.
29 Planar Point Location Sarnak and Tarjan. Sweep line + partially persistent binary search tree: Preprocessing time: O(n log n) Query time: O(log n) Space O(n) To get linear space: Balanced binary search tree with worst case O(1) memory modifications per update.
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